1. Field of the Invention
The present invention relates generally to a mobile communication system, and in particular, to an apparatus and method for coding/decoding block low density parity check (LDPC) codes.
2. Description of the Related Art
With the rapid development of mobile communication systems, it is necessary to develop technology capable of transmitting bulk data approximating the capacity of a wire network even in a wireless environment. To meet the increasing demand for a high-speed, high-capacity communication system capable of processing and transmitting various data such as image and radio data beyond the voice-oriented service, it is essential to increase the transmission efficiency of a system by using an appropriate channel coding scheme to thereby improve the overall system performance. However, the mobile communication system, because of its characteristics, inevitably generates errors during data transmission due to noise, interference and fading according to channel conditions. The generation of errors causes a loss of a great amount of information data.
In order to prevent the loss of the information data due to the generation of errors, various error control schemes are currently in use and are based in part on channel characteristics to thereby improve the reliability of the mobile communication system. The most typical error control scheme uses error correction codes.
With reference to FIG. 1, a description will now be made of a structure of a transmitter/receiver in a general mobile communication system.
FIG. 1 is a diagram illustrating a structure of a transmitter/receiver in a general mobile communication system. Referring to FIG. 1, a transmitter 100 includes an encoder 111, a modulator 113 and a radio frequency (RF) processor 115, and a receiver 150 includes an RF processor 151, a demodulator 153 and a decoder 155.
In the transmitter 100, transmission information data ‘u’, if generated, is delivered to the encoder 111. The encoder 111 generates a coded symbol ‘c’ by coding the information data ‘u’ with a predetermined coding scheme, and outputs the coded symbol ‘c’ to the modulator 113. The modulator 113 generates a modulation symbol ‘s’ by modulating the coded symbol ‘c’ with a predetermined modulation scheme, and outputs the modulation symbol ‘s’ to the RF processor 115. The RF processor 115 RF-processes the modulation symbol ‘s’ output from the modulator 113, and transmits the RF-processed signal over the air via an antenna ANT.
The signal transmitted over the air by the transmitter 100 in this way is received at the receiver 150 via its antenna ANT, and the signal received via the antenna is delivered to the RF processor 151. The RF processor 151 RF-processes the received signal, and outputs the RF-processed signal ‘r’ to the demodulator 153. The demodulator 153 demodulates the RF-processed signal ‘r’ output from the RF processor 151 using a demodulation scheme corresponding to the modulation scheme applied in the modulator 113, and outputs the demodulated signal ‘x’ to the decoder 155. The decoder 155 decodes the demodulated signal ‘x’ output from the demodulator 153 using a decoding scheme corresponding to the coding scheme applied in the encoder 111, and outputs the decoded signal ‘û’ as finally decoded information data.
In order for the receiver 150 to decode without errors the information data ‘u’ transmitted by the transmitter 100, there is a need for a high-performance encoder and decoder. Particularly, because a radio channel environment should be taken into consideration because of the characteristics of a mobile communication system, errors that can be generated due to the radio channel environment should be considered more seriously.
The most typical error correction codes include turbo codes and low density parity check (LDPC) codes.
It is well known that the turbo code is superior in performance gain to a convolutional code that is conventionally used for error correction, during high-speed data transmission. The turbo code is advantageous in that it can efficiently correct an error caused by noise generated in a transmission channel, thereby increasing the reliability of the data transmission. The LDPC code can be decoded using an iterative decoding algorithm base on a sum-product algorithm in a factor graph. Because a decoder for the LDPC code uses the sum-product algorithm-based iterative decoding algorithm, it is less complex than a decoder for the turbo code. In addition, the decoder for the LDPC code is easy to implement with a parallel processing decoder, compared with the decoder for the turbo code.
Shannon's channel coding theorem illustrates that reliable communication is possible only at a data rate not exceeding a channel capacity. However, Shannon's channel coding theorem has proposed no detailed channel coding/decoding method for supporting a data rate up to the maximum channel capacity limit. Generally, although a random code having a very large block size exhibits a performance approximating a channel capacity limit of Shannon's channel coding theorem, when a MAP (Maximum A Posteriori) or ML (Maximum Likelihood) decoding method is used, it is actually impossible to implement the decoding method because of its heavy calculation load.
The turbo code was proposed by Berrou, Glavieux and Thitimajshima in 1993, and exhibits a superior performance that approximates a channel capacity limit of Shannon's channel coding theorem. The proposal of the turbo code triggered active research on iterative decoding and graphical expression of codes, and LDPC codes proposed by Gallager in 1962 have been newly spotlighted in the research. Cycles exist in a factor graph of the turbo code and the LDPC code, and it is well known that iterative decoding in the factor graph of the LDPC code, where cycles exist, is suboptimal. Also, it has been experimentally proven that the LDPC code has excellent performance through iterative decoding. The LDPC code known to have the highest performance ever exhibits performances having a difference of only about 0.04 [dB] at a channel capacity limit of Shannon's channel coding theorem at a bit error rate (BER) 10−5, using a block size 107. In addition, although an LDPC code defined in Galois Field (GF) with q>2, i.e. GF(q), increases in complexity in its decoding process, it is much superior in performance to a binary code. However, there is no satisfactory theoretical description of successful decoding by an iterative decoding algorithm for the LDPC code defined in GF(q).
The LDPC code, proposed by Gallager, is defined by a parity check matrix in which most of elements have a value of 0 and minor elements except the elements having the value of 0 have a non-zero value, e.g., a value of 1. In the following description, it will be assumed that a non-zero value is a value of 1.
For example, an (N, j, k) LDPC code is a linear block code having a block length N, and is defined by a sparse parity check matrix in which each column has j elements having a value of 1, each row has k elements having a value of 1, and all of the elements except for the elements having the value of 1 have a value of 0.
An LDPC code in which a weight of each column in the parity check matrix is fixed to ‘j’ and a weight of each row in the parity check matrix is fixed to ‘k’ as described above, is called a “regular LDPC code.” Herein, the “weight” refers to the number of elements having a non-zero value among the elements constituting the parity check matrix. Unlike the regular LDPC code, an LDPC code in which the weight of each column in the parity check matrix and the weight of each row in the parity check matrix are not fixed is called an “irregular LDPC code.” It is generally known that the irregular LDPC code is superior in performance to the regular LDPC code. However, in the case of the irregular LDPC code, because the weight of each column and the weight of each row in the parity check matrix are not fixed, i.e. are irregular, the weight of each column in the parity check matrix and the weight of each row in the parity check matrix must be properly adjusted in order to guarantee the superior performance.
With reference to FIG. 2, a description will now be made of a parity check matrix of an (8, 2, 4) LDPC code as an example of an (N, j, k) LDPC code.
FIG. 2 is a diagram illustrating a parity check matrix of a general (8, 2, 4) LDPC code. Referring to FIG. 2, a parity check matrix H of the (8, 2, 4) LDPC code is comprised of 8 columns and 4 rows, wherein a weight of each column is fixed to 2 and a weight of each row is fixed to 4. Because the weight of each column and the weight of each row in the parity check matrix are regular as stated above, the (8, 2, 4) LDPC code illustrated in FIG. 2 becomes a regular LDPC code.
The parity check matrix of the (8, 2, 4) LDPC code has been described so far with reference to FIG. 2. Next, a factor graph of the (8, 2, 4) LDPC code described in connection with FIG. 2 will be described herein below with reference to FIG. 3.
FIG. 3 is a diagram illustrating a factor graph of the (8, 2, 4) LDPC code of FIG. 2. Referring to FIG. 3, a factor graph of the (8, 2, 4) LDPC code is comprised of 8 variable nodes of x1 300, x2 302, x3 304, x4 306, x5 308, x6 310, x7 312 and x8 314, and 4 check nodes 316, 318, 320 and 322. When an element having a value of 1, i.e. a non-zero value, exists at a point where an ith row and a jth column of the parity check matrix of the (8, 2, 4) LDPC code cross each other, a branch is created between a variable node xi and a jth check node.
Because the parity check matrix of the LDPC code has a very small weight as described above, it is possible to perform decoding through iterative decoding even in a block code having a relatively long length, that exhibits performance approximating a channel capacity limit of Shannon's channel coding theorem, such as a turbo code, while continuously increasing a block length of the block code. MacKay and Neal have proven that an iterative decoding process of an LDPC code using a flow transfer scheme approximates an iterative decoding process of a turbo code in performance.
In order to generate a high-performance LDPC code, the following conditions should be satisfied.
(1) Cycles on a factor graph of an LDPC code should be considered.
The term “cycle” refers to a loop formed by the edges connecting the variable nodes to the check nodes in a factor graph of an LDPC code, and a length of the cycle is defined as the number of edges constituting the loop. A long cycle means that the number of edges connecting the variable nodes to the check nodes constituting the loop in the factor graph of the LDPC code is large. In contrast, a short cycle means that the number of edges connecting the variable nodes to the check nodes constituting the loop in the factor graph of the LDPC code is small.
As cycles in the factor graph of the LDPC code become longer, the performance efficiency of the LDPC code increases, for the following reasons. That is, when long cycles are generated in the factor graph of the LDPC code, it is possible to prevent performance degradation such as an error floor occurring when too many cycles with a short length exist on the factor graph of the LDPC code.
(2) Efficient coding of an LDPC code should be considered.
It is difficult for the LDPC code to undergo real-time coding compared with a convolutional code or a turbo code because of its high coding complexity. In order to reduce the coding complexity of the LDPC code, a Repeat Accumulate (RA) code has been proposed. However, the RA code also has a limitation in reducing the coding complexity of the LDPC code. Therefore, efficient coding of the LDPC code should be taken into consideration.
(3) Degree distribution on a factor graph of an LDPC code should be considered.
Generally, an irregular LDPC code is superior in performance to a regular LDPC code, because a factor graph of the irregular LDPC code has various degrees. The term “degree” refers to the number of edges connected to the variable nodes and the check nodes in the factor graph of the LDPC code. Further, the phrase “degree distribution” on a factor graph of an LDPC code refers to a ratio of the number of nodes having a particular degree to the total number of nodes. It has been proven by Richardson that an LDPC code having a particular degree distribution is superior in performance.
With reference to FIG. 4, a description will now be made of a parity check matrix of a block LDPC code.
FIG. 4 is a diagram illustrating a parity check matrix of a general block LDPC code. Before a description of FIG. 4 is given, it should be noted that the block LDPC code is a new LDPC code for which not only efficient coding but also efficient storage and performance improvement of a parity check matrix were considered, and the block LDPC code is an LDPC code extended by generalizing a structure of a regular LDPC code. Referring to FIG. 4, a parity check matrix of the block LDPC code is divided into a plurality of partial blocks, and a permutation matrix is mapped to each of the partial blocks. In FIG. 4, ‘P’ represents a permutation matrix having an Ns×Ns size, and a superscript (or exponent) apq of the permutation matrix P is either 0≦apq≦Ns−1 or apq=∞.
In addition, ‘p’ indicates that a corresponding permutation matrix is located in the pth row of the partial blocks of the parity check matrix, and ‘q’ indicates that a corresponding permutation matrix is located in the qth column of the partial blocks of the parity check matrix. That is, Pαpq represents a permutation matrix located in a partial block where the pth row and the qth column of the parity check matrix comprised of a plurality of partial blocks cross each other. That is, the ‘p’ and the ‘q’ represent the number of rows and the number of columns of partial blocks corresponding to an information part in the parity check matrix, respectively.
The permutation matrix will now be described with reference to FIG. 5.
FIG. 5 is a diagram illustrating the permutation matrix P of FIG. 4. As illustrated in FIG. 5, the permutation matrix P is a square matrix having an Ns×Ns, size, and each of Ns columns constituting the permutation matrix P has a weight of 1 and each of Ns rows constituting the permutation matrix P also has a weight of 1. Herein, although a size of the permutation matrix P is expressed as Ns×Ns, it will also be expressed as Ns because the permutation matrix P is a square matrix.
In FIG. 4, a permutation matrix P with a superscript apq=0, i.e. a permutation matrix P0, represents an identity matrix INs×Ns, and a permutation matrix P with a superscript apq=∞, i.e. a permutation matrix P∞, represents a zero matrix. Herein, INs×Ns represents an identity matrix with a size Ns×Ns.
In the entire parity check matrix of the block LDPC code illustrated in FIG. 4, because the total number of rows is Ns×p and the total number of columns is Ns×q(for p≦q), when the entire parity check matrix of the LDPC code has a full rank, a coding rate can be expressed as Equation (1) regardless of a size of the partial blocks.
                    R        =                                                                              N                  s                                ×                q                            -                                                N                  s                                ×                p                                                                    N                s                            ×              q                                =                                                    q                -                p                            q                        =                          1              -                              p                q                                                                        (        1        )            
If apq≠∞ for all p and q, the permutation matrixes corresponding to the partial blocks are not zero matrixes, and the partial blocks constitute a regular LDPC code in which the weight value of each column and the weight value of each row in each of the permutation matrixes corresponding to the partial blocks are p and q, respectively. Herein, each of permutation matrixes corresponding to the partial blocks will be referred to as a “partial matrix.”
Because (p-1) dependent rows exist in the entire parity check matrix, a coding rate is greater than the coding rate calculated by Equation (1). In the case of the block LDPC code, if a weight position of a first row of each of the partial matrixes constituting the entire parity check matrix is determined, the weight positions of the remaining (Ns-1) rows can be determined. Therefore, the required size of a memory is reduced to 1/Ns as compared with the case where the weights are irregularly selected to store information on the entire parity check matrix.
As described above, the term “cycle” refers to a loop formed by the edges connecting the variable nodes to the check nodes in a factor graph of an LDPC code, and a length of the cycle is defined as the number of edges constituting the loop. A long cycle means that the number of edges connecting the variable nodes to the check nodes constituting the loop in the factor graph of the LDPC code is large. As cycles in the factor graph of the LDPC code become longer, the performance efficiency of the LDPC code increases.
In contrast, as cycles in the factor graph of the LDPC code become shorter, an error correction capability of the LDPC code decreases because performance degradation such as an error floor occurs. That is, when there are many cycles with a short length in a factor graph of the LDPC code, information on a particular node belonging to the cycle with a short length, starting therefrom, returns after a small number of iterations. As the number of iterations increases, the information returns to the corresponding node more frequently, so that the information cannot be correctly updated, thereby causing a deterioration in an error correction capability of the LDPC code.
With reference to FIG. 6, a description will now be made of a cycle structure of a block LDPC code.
FIG. 6 is a diagram illustrating a cycle structure of a block LDPC code of which a parity check matrix is comprised of 4 partial matrixes. Before a description of FIG. 6 is given, it should be noted that the block LDPC code is a new LDPC code for which not only efficient coding but also efficient storage and performance improvement of a parity check matrix were considered. The block LDPC code is also an LDPC code extended by generalizing a structure of a regular LDPC code. A parity check matrix of the block LDPC code illustrated in FIG. 6 is comprised of 4 partial blocks, a diagonal line represents a position where the elements having a value of 1 are located, and the portions other than the diagonal-lined portions represent positions where the elements having a value of 0 are located. In addition, ‘P’ represents the same permutation matrix as the permutation matrix described in conjunction with FIG. 5.
In order to analyze a cycle structure of the block LDPC code illustrated in FIG. 6, an element having a value of 1 located in an ith row of a partial matrix pa is defined as a reference element, and an element having a value of 1 located in the ith row will be referred to as a “0-point.” Herein, “partial matrix” will refer to a matrix corresponding to the partial block. The 0-point is located in an (i+a)th column of the partial matrix Pa.
An element having a value of 1 in a partial matrix Pb, located in the same row as the 0-point, will be referred to as a “1-point.” For the same reason as the 0-point, the 1-point is located in an (i+b)th column of the partial matrix Pb.
Next, an element having a value of 1 in a partial matrix Pc, located in the same column as the 1-point, will be referred to as a “2-point.” Because the partial matrix Pc is a matrix acquired by shifting respective columns of an identity matrix I to the right with respect to a modulo Ns by c, the 2-point is located in an (i+b−c)th row of the partial matrix Pc.
In addition, an element having a value of 1 in a partial matrix Pd, located in the same row as the 2-point, will be referred to as a “3-point.” The 3-point is located in an (i+b−c+d)th column of the partial matrix Pd.
Finally, an element having a value of 1 in the partial matrix Pa, located in the same column as the 3-point, will be referred to as a “4-point.” The 4-point is located in an (i+b−c+d−a)th row of the partial matrix Pa.
In the cycle structure of the LDPC code illustrated in FIG. 6, if a cycle with a length of 4 exists, the 0-point and the 4-point are located in the same position. That is, a relation between the 0-point and the 4-point is defined by Equation (2)i≅i+b−c+d−a(mod Ns) ori+a≅i+b−c+d(mod Ns)  (2)
Equation (2) can be rewritten as Equation (3)a+c≅b+d(mod Ns)  (3)
As a result, when the relationship of Equation (3) is satisfied, a cycle with a length 4 is generated. Generally, when a 0-point and a 4p-point are first identical to each other, a relation of i≅i+p(b−c+d−e)(mod Ns) is given, and the following relation shown in Equation (4) is satisfied.p(a−b+c−d)≅0(mod Ns)  (4)
In other words, if a positive integer having a minimum value among the positive integers satisfying Equation (4) for a given a, b, c and d is defined as ‘p’, a cycle with a length of 4p becomes a cycle having a minimum length in the cycle structure of the block LDPC code illustrated in FIG. 6.
In conclusion, as described above, for (a−b+c−d)≠0, if gcd(Ns, a−b+c−d)=1 is satisfied, then p=Ns. Herein, the gcd(Ns, a−b+c−d) is the function for calculating the “greatest common divisor” of the integers Ns and a−b+c−d. Therefore, a cycle with a length of 4Ns becomes a cycle with a minimum length.
A Richardson-Urbanke technique will be used as a coding technique for the block LDPC code. Because the Richardson-Urbanke technique is used as a coding technique, coding complexity can be minimized as the form of a parity check matrix becomes similar to the form of a full lower triangular matrix.
With reference to FIG. 7, a description will now be made of a parity check matrix having a form similar to the form of the full lower triangular matrix.
FIG. 7 is a diagram illustrating a parity check matrix having a form similar to the form of the full lower triangular matrix. The parity check matrix illustrated in FIG. 7 is different from the parity check matrix having a form of the full lower triangular matrix in the form of the parity part. In FIG. 7, a superscript (or exponent) apq of the permutation matrix P of an information part is either 0≦apq≦Ns−1 or apq=∞, as described above. A permutation matrix P with a superscript apq=0, i.e. a permutation matrix P0, of the information part represents an identity matrix INs×Nx, and a permutation matrix P with a superscript apq∞, i.e. a permutation matrix P∞, represents a zero matrix. In. FIG. 7, ‘p’ represents the number of rows of partial blocks mapped to the information part, and ‘q’ represents the number of columns of partial blocks mapped to the parity part. Also, superscripts ap, x and y of the permutation matrixes P mapped to the parity part represent exponents of the permutation matrix P. However, for the convenience of explanation, the different superscripts ap, x and y are used to distinguish the parity part from the information part. That is, in FIG. 7, pal and Pap are also permutation matrixes, and the superscripts a1 to ap are sequentially indexed to partial matrixes located in a diagonal part of the parity part. In addition, Px and Py are also permutation matrixes, and for the convenience of explanation, they are indexed in a different way to distinguish the parity part from the information part. If a block length of a block LDPC code having the parity check matrix illustrated in FIG. 7 is assumed to be N, the coding complexity of the block LDPC code linearly increases with respect to the block length N (0(N)).
The biggest problem of the LDPC code having the parity check matrix of FIG. 7 is that if a length of a partial block is defined as Ns, Ns check nodes whose degrees are always 1 in a factor graph of the block LDPC code are generated. The check nodes with a degree of 1 cannot affect the performance improvement based on the iterative decoding. Therefore, a standard irregular LDPC code based on the Richardson-Urbanke technique does not include a check node with a degree of 1. Therefore, a parity check matrix of FIG. 7 will be assumed as a basic parity check matrix in order to design a parity check matrix such that it enables efficient coding while not including a check node with a degree of 1. In the parity check matrix of FIG. 7 comprised of the partial matrixes, the selection of a partial matrix is a very important factor for a performance improvement of the block LDPC code, so that finding an appropriate selection criterion for the partial matrix also becomes a very important factor.
A description will now be made of a method for designing the parity check matrix of the block LDPC code based on the foregoing block LDPC code.
In order to facilitate a method of designing a parity check matrix of the block LDPC code and a method for coding the block LDPC code, the parity check matrix illustrated in FIG. 7 is assumed to be formed with 6 partial matrixes as illustrated in FIG. 8.
FIG. 8 is a diagram illustrating the parity check matrix of FIG. 7 which is divided into 6 partial blocks. Referring to FIG. 8, a parity check matrix of the block LDPC code illustrated in FIG. 7 is divided into an information part ‘s’, a first parity part p1, and a second parity part p2. The information part ‘s’ represents a part of the parity check matrix, mapped to an actual information word during the process of coding a block LDPC code, like the information part described in conjunction with FIG. 7, but for the convenience of explanation, the information part ‘s’ is represented by different reference letters. The first parity part p1 and the second parity part p2 represent a part of the parity check matrix, mapped to an actual parity during the process of coding the block LDPC code, like the parity part described in conjunction with FIG. 7, and the parity part is divided into two parts.
Partial matrixes A and C correspond to partial blocks A (802) and C (804) of the information part ‘s’, partial matrixes B and D correspond to partial blocks B (806) and D (808) of the first parity part p1, and partial matrixes T and E correspond to partial blocks T (810) and E (812) of the second parity part p2. Although the parity check matrix is divided into 7 partial blocks in FIG. 8, it should be noted that ‘0’ is not a separate partial block and because the partial matrix T corresponding to the partial block T (810) have a full lower triangular form, a region where zero matrixes are arranged on the basis of a diagonal is represented by ‘0’. A process of simplifying a coding method using the partial matrixes of the information part ‘s’, the first parity part p1 and the second parity part p2 will be described later with reference to FIG. 10.
The partial matrixes of FIG. 8 will now be described herein below with reference to FIG. 9.
FIG. 9 is a diagram illustrating a transpose matrix of the partial matrix B shown in FIG. 8, the partial matrix E, the partial matrix T, and an inverse matrix of the partial matrix T, in the parity check matrix of FIG. 7.
Referring to FIG. 9, a partial matrix BT represents a transpose matrix of the partial matrix B, and a partial matrix T−1 represents an inverse matrix of the partial matrix T. The P(k1˜k2) represents
            ∏              i        =                  k          i                            k        2              ⁢          P              a        i              =            P                        ∑                      i            =                          k              i                                            k            2                          ⁢                  a          i                      .  The permutation matrixes illustrated in FIG. 9, for example, Pa1, may be an identity matrix. As described above, if a superscript of the permutation matrix, i.e. a1 is 0, the Pa1 will be an identity matrix. Also, if a superscript of the permutation matrix, i.e. a, increases by a predetermined value, the permutation matrix is cyclic shifted by the predetermined value, so the permutation matrix Pa1 will be an identity matrix.
With reference to FIG. 10, a description will now be made of a process of designing a parity check matrix of the block LDPC code.
FIG. 10 is a flowchart illustrating a procedure for generating a parity check matrix of a general block LDPC code. Before a description of FIG. 10 is given, it should be noted that in order to generate a block LDPC code, a codeword size and a coding rate of a block LDPC code to be generated must be determined, and a size of a parity check matrix must be determined according to the determined codeword size and coding rate. If a codeword size of the block LDPC code is represented by N and a coding rate is represented by R, a size of a parity check matrix becomes N(1−R)×N. Actually, the procedure for generating a parity check matrix of a block LDPC code illustrated in FIG. 10 is performed only once, because the parity check matrix is initially generated to be suitable for a situation of a communication system and thereafter, the generated parity check matrix is used.
Referring to FIG. 10, in step 1011, a controller divides a parity check matrix with the size N(1−R)×N into a total of p×q blocks, including p blocks in a horizontal axis and q blocks in a vertical axis, and then proceeds to step 1013. Because each of the blocks has a size of Ns×Ns, the parity check matrix is comprised of Ns×p columns and Ns×q rows. In step 1013, the controller classifies the p×q blocks divided from the parity check matrix into an information part ‘s’, a first parity part p1, and a second parity part p2, and then proceeds to steps 1015 and 1021.
In step 1015, the controller separates the information part ‘s’ into non-zero blocks, or non-zero matrixes, and zero blocks, or zero matrixes according to the degree of distribution for guaranteeing a good performance of the block LDPC code, and then proceeds to step 1017. Because the degree of distribution for guaranteeing a good performance of the block LDPC code has been described above, a detailed description thereof will omitted herein. In step 1017, the controller determines the permutation matrixes Papq such that a minimum cycle length of a block cycle should be maximized as described above in the non-zero matrix portions in blocks having a low degree from among the blocks determined according to the degree distribution for guaranteeing a good performance of the block LDPC code, and then proceeds to step 1019. The permutation matrixes Papq should be determined taking into consideration the block cycles of not only the information part ‘s’ but also the first parity part p1 and the second parity part p2.
In step 1019, the controller randomly determines the permutation matrixes Papq in the non-zero matrix portions in the blocks having a high degree among the blocks determined according to the degree distribution for guaranteeing a good performance of the block LDPC code, and then ends the procedure. Even when the permutation matrixes Papq to be applied to the non-zero matrix portions in the blocks having a high degree are determined, the permutation matrixes Papq must be determined by such that a minimum cycle length of a block cycle is maximized, and the permutation matrixes Papq are determined considering the block cycles of not only the information part ‘s’ but also the first parity part p1 and the second parity part p2. An example of the permutation matrixes Papq arranged in the information part ‘s’ of the parity check matrix is illustrated in FIG. 7.
In step 1021, the controller divides the first part p1 and the second parity part p2 into 4 partial matrixes B, T, D and E, and then proceeds to step 1023. In step 1023, the controller inputs the non-zero permutation matrixes PY and Pa1 to 2 partial blocks among the partial blocks constituting the partial matrix B, and then proceeds to step 1025. The structure for inputting the non-zero permutation matrixes Py and Pa1 to 2 partial blocks among the partial blocks constituting the partial matrix B has been described with reference to FIG. 9.
In step 1025, the controller inputs the identity matrixes I to the diagonal partial blocks of the partial matrix T, inputs the particular permutation matrixes Pa2, Pa3, . . . , Pam-1 to (i, i+1)th partial blocks under the diagonal components of the partial matrix T, and then proceeds to step 1027. The structure for inputting the identity matrixes I to the diagonal partial blocks of the partial matrix T and inputting the particular permutation matrixes Pa2, Pa3, . . . , Pam-1 to (i, i+1)th partial blocks under the diagonal components of the partial matrix T has been described with reference to FIG. 9.
In step 1027, the controller inputs a partial matrix Px to the partial matrix D, and then proceeds to step 1029. In step 1029, the controller inputs a permutation matrix Pam to only the last partial block in the partial matrix E, and then ends the procedure. The structure for inputting the 2 permutation matrixes Pam to only the last partial block among the partial blocks constituting the partial matrix E has been described with reference to FIG. 9.